Operations Research

Motivation

Set covering asks for the fewest members of a family of sets whose union is everything the family covers. It models crew scheduling, facility siting, test-suite reduction, logic minimization and fault testing; Johnson names the last two as its practical applications. Karp showed in 1972 that the decision version is NP-complete (Karp 1972), so in practice one runs a heuristic and asks how far from optimal it can be.

David S. Johnson's 1974 paper Approximation Algorithms for Combinatorial Problems (JCSS 9, 256–278) is one of the founding papers of the worst-case analysis of approximation algorithms. For set covering it analyses the obvious greedy rule, repeatedly take a set that covers the most still-uncovered points, and proves that on families whose sets have at most $k$ elements its output is never more than the harmonic number $H(k) = \sum_{j=1}^k 1/j$ times the optimum, and that this factor is attained.

Timeline.

• 1974: Johnson proves the $H(k)$ bound for unweighted set cover with sets of size at most $k$, together with a matching family of examples (this mission).
• 1975: Lovász proves the same bound for the fractional relaxation, giving an integrality-gap statement (Lovász 1975).
• 1979: Chvátal extends the bound to weighted set cover, with the greedy rule choosing the set of least cost per newly covered point (Chvátal 1979).
• 1998: Feige shows that no polynomial-time algorithm achieves $(1-\varepsilon)\ln n$ unless NP has slightly superpolynomial deterministic algorithms (Feige 1998), so the greedy guarantee is essentially the best possible.

Setting

An input $F$ of SET COVERING I is a finite family $\{S_1, \dots, S_p\}$ of finite sets. The set to be covered is $T = \bigcup_{S \in F} S$. A subcover is a subfamily $F' \subseteq F$ with $\bigcup_{S \in F'} S = T$, and its measure is $|F'|$. The optimum $F^*$ is the minimum measure of a subcover; $F$ itself is a subcover, so the minimum exists. The subproblem SC(k) restricts the inputs to families no set of which has more than $k$ elements.

Algorithm C1 keeps a family SUB of chosen sets, the set UNCOV of uncovered points, and an array SET$[i]$ holding the still-uncovered part of $S_i$. It starts with SUB $= \emptyset$, UNCOV $= T$, SET$[i] = S_i$. While UNCOV is nonempty it chooses an index $j$ with $|\mathrm{SET}[j]|$ maximal, adds $S_j$ to SUB, and removes SET$[j]$ from UNCOV and from every SET$[i]$. When UNCOV is empty it returns SUB. When several indices tie at Step 3 any of them may be chosen, so one input can have several choosable outputs. Following Section 2 of the paper, the algorithm's value $C1(F)$ is the worst choosable output, here the largest, and the ratio is $r(C1, F) = C1(F)/F^*$.

For the proof the paper introduces configurations $K = \langle N_K, \mathrm{UNCOV}_K, \langle \mathrm{SET}_K[1], \dots, \mathrm{SET}_K[N_K]\rangle\rangle$ with $\bigcup_i \mathrm{SET}_K[i] = \mathrm{UNCOV}_K$, runs from a configuration (sequences of admissible choices ending when UNCOV is empty), $\mathrm{Numbers}(R)$, the set of indices chosen in a run $R$, and calls a set $M$ selectable from $K$ if $M = \mathrm{Numbers}(R)$ for some run $R$ from $K$. Write $n(K, i) = |\mathrm{SET}_K[i]|$.

Formalization targets

Goal: Theorem 4

For every $k \ge 1$:

$$\text{for every input } F \in SC(k) \text{ and every choosable } F_1:\quad |F_1| \le H(k)\cdot F^*,$$ $$\text{and some } F \in SC(k) \text{ with } F^* > 0 \text{ has a choosable } F_1 \text{ with } |F_1| = H(k)\cdot F^*.$$

The paper states this as $R[C1, SC(k)](n) \le \sum_{j=1}^k (1/j)$ for all $n > 0$, with equality for all sufficiently large $n$. The two-part form above is the size-free equivalent.

Milestones

1. Lemma 1. For a subcover $F_1$ with index set $M1 = \{i : S_i \in F_1\}$ and $K$ the configuration after Step 1: $F_1$ is choosable by C1 if and only if $M1$ is selectable from $K$.
2. Lemma 2. For any configuration $K$, any $M1$ selectable from $K$ and any $M0$ with $\bigcup_{i \in M0} \mathrm{SET}_K[i] = \mathrm{UNCOV}_K$:

$$|M1| \le \sum_{i \in M0} \sum_{j=1}^{n(K,i)} \frac{1}{j}.$$

3. Fig. 1. For every $k \ge 1$ there is an explicit input of $SC(k)$ on $k \cdot k!$ points with $F^* = k!$ and a choosable output of $k!\,H(k)$ sets.

Significance

The result. Theorem 4 is the first proof that greedy set cover has a worst-case guarantee depending only on the largest set size, and it pins the guarantee down exactly: the constant $H(k)$ cannot be lowered for any $k$. Since $H(k) \le 1 + \ln k$, it also gives the well-known $1 + \ln n$ bound for general inputs. The $H(k)$ bound and its later refinements are the standard reference point for analyses of greedy covering, dual fitting and submodular covering.

Formalizing it. The theorem has been proved since 1974. As far as a search of the platform shows, no machine-checked proof of it exists: the platform holds a Kearns–Vazirani-style statement ComputationalLearning.greedy_set_cover (the $\mathrm{opt}\cdot\ln|U|$ form for a greedy sequence, still open) and a dual-fitting certificate lemma for weighted set cover, neither of which covers the $SC(k)$ bound, the tie-breaking semantics or the tightness construction. A complete development provides both halves of Theorem 4, the configuration and run machinery of Lemmas 1–2, and the explicit Fig. 1 family.

Difficulty

The obvious argument charges each chosen set to the points it newly covers and compares the charges with an optimal cover. A statement about the initial input alone, with the original sizes of the optimal sets, does not survive a single greedy step: after a step the optimal sets are only partly uncovered and the remaining run faces a different instance. This is why Lemma 2 is stated for an arbitrary configuration, in terms of the current sizes $n(K, i)$, and for an arbitrary covering subfamily $M0$. Because Step 3 breaks ties arbitrarily, the statement must hold for every admissible run, and a formalization that fixes one tie-breaking rule proves a weaker upper bound and cannot express the tightness example, which relies on adversarial ties at every stage.

For the tightness half, the difficulty is bookkeeping: showing that the $k!/j$ blocks of each segment are admissible choices at each stage and that no cover uses fewer than $k!$ sets.

Formalization scope

• An input is an indexed family S : ι → Finset α over a finite index type ι and a ground type with decidable equality. The indices play the role of $1, \dots, N$; two indices may carry the same set, which only widens the input class. The family, subcovers and $F^*$ are taken over the set of sets family S, as on the page. $F^*$ is a Finset.inf' over the nonempty finite set of subcovers; if $T = \emptyset$ then $F^* = 0$.
• C1 is a nondeterministic step relation: a step is allowed for every index maximizing $|\mathrm{SET}[j]|$. An output is choosable if a finite chain of steps from the initial state reaches a halting state with that SUB. No tie-breaking rule is fixed.
• The paper's $R[A, P](n)$ is a maximum over inputs of size at most $n$ in an unspecified notation; it is replaced by the size-free two-part statement above, which is equivalent because $R$ is a maximum over finitely many inputs and nondecreasing in $n$.
• Ratios are stated multiplicatively in $\mathbb{Q}$ ($|F_1| \le H(k)\cdot F^*$), never as a quotient, so an input with $F^* = 0$ does not make the bound vacuous, and the attainment part requires $F^* > 0$. $H(k)$ is Mathlib's harmonic k.
• Configurations carry the covering condition as a field; runs are an inductive predicate on the list of chosen indices; Selectable K M means $M$ is the set of indices of some run.
• Lemma 1 assumes the family's sets are pairwise distinct (the paper's family is a set of sets); without that the index set $\{i : S_i \in F_1\}$ may contain a duplicate index C1 never chose.
• Trivializing formalizations are ruled out: a deterministic tie-break, a ratio written as a division, the original set sizes in place of $n(K, i)$ in Lemma 2, or an attaining input with $F^* = 0$ would each change the theorem.

Contributions welcome: proofs of Lemma 2 (the core induction), of Lemma 1, of the Fig. 1 run, and of Theorem 4 from these; the configuration/run layer and the Fig. 1 family are reusable for other greedy covering analyses.

Selected references

• David S. Johnson, Approximation algorithms for combinatorial problems, Journal of Computer and System Sciences 9 (1974), 256–278. https://doi.org/10.1016/S0022-0000(74)80044-9
• Richard M. Karp, Reducibility among combinatorial problems, in Complexity of Computer Computations, Plenum, 1972, 85–103. https://doi.org/10.1007/978-1-4684-2001-2_9
• László Lovász, On the ratio of optimal integral and fractional covers, Discrete Mathematics 13 (1975), 383–390. https://doi.org/10.1016/0012-365X(75)90058-8
• Vašek Chvátal, A greedy heuristic for the set-covering problem, Mathematics of Operations Research 4 (1979), 233–235. https://doi.org/10.1287/moor.4.3.233
• Uriel Feige, A threshold of ln n for approximating set cover, Journal of the ACM 45 (1998), 634–652. https://doi.org/10.1145/285055.285059

Motivation

Maximum satisfiability asks for a truth assignment satisfying as many clauses of a propositional formula as possible. The paper notes that the restriction $MS(k)$, in which every clause has at least $k$ literals, is polynomial complete for every $k \ge 1$, so exact optimization is out of reach in general and one asks instead how close a fast algorithm is guaranteed to come. David S. Johnson's 1974 paper Approximation Algorithms for Combinatorial Problems (J. Comput. System Sci. 9, 256–278) set up a framework for exactly this question — optimization problems, nondeterministic approximation algorithms, and the worst-case ratio between the optimum and the algorithm's output — and applied it to subset-sum, maximum satisfiability, set covering, graph coloring and maximum clique. It is one of the founding papers of the theory of approximation algorithms.

Section 4 of the paper treats maximum satisfiability with two algorithms. This mission covers the first, a greedy literal-selection rule called B1, and its exact worst-case ratio (Theorem 2). A companion mission covers the weighted algorithm B2 (Theorem 3).

Timeline, for orientation:

• 1971–1972: Cook and Karp establish NP-completeness of satisfiability and of many combinatorial problems.
• 1974: Johnson proves that B1 has worst-case ratio exactly $(k+1)/k$ on $MS(k)$ and that the weighted algorithm B2 achieves $2^k/(2^k-1)$ (Theorems 2 and 3).
• 1990s: semidefinite and LP-based algorithms (Goemans–Williamson, SIAM J. Discrete Math. 1994) improve the constants for general MAX-SAT.

Setting

Let $L = \bigcup_{i>0}\{x_i, \bar x_i\}$ be the set of literals; the complement of $x_i$ is $\bar x_i$ and conversely. A clause is a finite set $C \subseteq L$. A truth assignment is a set $T \subseteq L$ containing no complementary pair $\{x_i, \bar x_i\}$; it may leave variables unassigned. $T$ satisfies $C$ if $C \cap T \ne \emptyset$.

An input is a finite set $S$ of clauses. Its feasible solutions are the subsets $S' \subseteq S$ satisfied by a single truth assignment, measured by $|S'|$, and the optimum is

$$S^* = \max\{|S'| : S' \subseteq S,\ \text{some truth assignment satisfies every } C \in S'\}.$$

The subproblem $MS(k)$ admits only inputs whose clauses each contain at least $k$ distinct literals.

Algorithm B1 keeps four variables: SUB (clauses already satisfied), LEFT (clauses not yet satisfied), TRUE (literals made true) and LIT (literals still available). It starts with SUB $=$ TRUE $= \emptyset$, LEFT $= S$, LIT $= L$. While some literal of LIT occurs in a clause of LEFT, it picks a literal $y \in$ LIT contained in the most clauses of LEFT, moves those clauses $YT$ from LEFT to SUB, adds $y$ to TRUE, and removes $y$ and $\bar y$ from LIT. When no literal of LIT occurs in LEFT it returns SUB.

The choice of $y$ is not determined when several literals tie. Following the paper's framework, every output reachable by some sequence of admissible choices is choosable, and the performance of B1 on $S$ is the smallest $|X|$ over choosable outputs $X$. The worst-case ratio on inputs of size at most $n$ is

$$R[B1, MS(k)](n) = \max\{S^*/B1(S) : S \in MS(k),\ |S| \le n\}.$$

Formalization targets

Goal: Theorem 2 (p. 262)

For all $k \ge 1$,

$$R[B1, MS(k)](n) \le \frac{k+1}{k}\quad\text{for all } n > 0,$$

with equality for all sufficiently large $n$. In the size-free form used here: every choosable output $X$ on every $S \in MS(k)$ satisfies $k\,S^* \le (k+1)\,|X|$, and for every $k \ge 1$ some $S \in MS(k)$ has a choosable $X$ with $|X| > 0$ and $k\,S^* = (k+1)\,|X|$.

Milestones (from the proof of Theorem 2, pp. 262–263)

1. In each iteration, the number of clauses saved (added to SUB) is at least the number of clauses remaining in LEFT that are wounded (lose a literal from LIT without being satisfied).
2. When B1 halts, every clause left in LEFT is dead: each of its literals has had its complement made true.
3. When B1 halts on an input of $MS(k)$, $|\mathrm{SUB}| \ge k\,|\mathrm{LEFT}|$, and SUB and LEFT partition $S$.
4. On the four-clause input $\{\{x_1,x_2,x_3\},\{\bar x_1,x_4,x_5\},\{\bar x_2,x_6,x_7\},\{\bar x_3,x_8,x_9\}\}$ of $MS(3)$, $S^* = 4$ while B1 may return three clauses.

Significance

The bound is stronger than a ratio: milestone 3 shows that B1 always satisfies at least $\tfrac{k}{k+1}|S|$ clauses, whatever the optimum. The tightness half shows that this simple greedy rule cannot be analysed any better, which is what motivated the weighted algorithm B2 of the same section, with ratio $2^k/(2^k-1)$. The pair of theorems is an early instance of a now standard pattern: a potential-style counting argument for an upper bound, and an adversarial tie-breaking instance for the matching lower bound.

The result is proved in the paper; it has not, to our knowledge, been machine-checked. This mission produces a formal model of Johnson's framework for a maximization problem with a nondeterministic algorithm, a formal proof of the upper bound through the "saved versus wounded" accounting, and explicit tightness instances for every $k \ge 1$. The paper spells out only $k = 3$ and states that "similar examples can be constructed for any other $k > 0$"; the formal goal requires them for all $k$.

Difficulty

The upper bound needs an invariant over entire runs, not over a single step: a clause wounded in one iteration may be saved in a later one, so wounds and saves must be tallied globally, and the count of wounds received by a clause that ends in LEFT must be matched with its number of literals. That matching relies on the facts that B1 never makes both a literal and its complement true and that a clause containing a true literal has already left LEFT. Clauses containing both $x_i$ and $\bar x_i$ are allowed and have to be handled.

The lower bound cannot be obtained from a fixed tie-breaking rule: the attaining run chooses negative literals whose count merely ties the maximum. For general $k$ the instance has to be built so that every literal occurs in few enough clauses that the adversarial choice is admissible at every step; at $k = 1$ the paper's pattern degenerates and needs adjusting.

Formalization scope

• A literal is a pair (variable index in $\mathbb N$, sign); a clause is a Finset of literals; an input is a Finset of clauses, so duplicate clauses are not allowed, as on the page. Tautological clauses are allowed.
• A truth assignment is a Set of literals without a complementary pair (partial, as in the paper). $S^*$ is the maximum of $|S'|$ over the finite nonempty family of satisfiable subsets, taken with Finset.sup'.
• B1 is a nondeterministic run relation: a state holds SUB, LEFT, TRUE and the set of decided variables (LIT is its complement, since $L$ is infinite); one step chooses any literal of LIT, of either sign, with maximum count; "choosable" is reachability of a halting state with the given SUB. No tie-break is fixed. A formalization that picks a variable and then its better sign, or that resolves ties deterministically, is a different algorithm and would make the tightness half false.
• The ratio $R[B1, MS(k)](n)$, whose problem size is left unspecified in the paper, is replaced by its size-free equivalent, and ratios are written multiplicatively in $\mathbb N$: $k\,S^* \le (k+1)|X|$. The tightness half requires $|X| > 0$, so the empty input cannot witness it.
• The running time $O(n \log n)$ is not stated.

Welcome contributions: proofs of the milestones, the invariants of reachable B1 states (SUB and LEFT partition $S$; TRUE is consistent and exactly covers the decided variables; no clause of LEFT meets TRUE), and the family of tightness instances for general $k$. The run-relation encoding of choosable outputs is reusable for the other algorithms of the paper.

Selected references

• D. S. Johnson, Approximation algorithms for combinatorial problems, Journal of Computer and System Sciences 9 (1974), 256–278. https://doi.org/10.1016/S0022-0000(74)80044-9
• R. M. Karp, Reducibility among combinatorial problems, in Complexity of Computer Computations, Plenum, 1972, 85–103. https://doi.org/10.1007/978-1-4684-2001-2_9
• M. X. Goemans and D. P. Williamson, New 3/4-approximation algorithms for the maximum satisfiability problem, SIAM Journal on Discrete Mathematics 7 (1994), 656–666. https://doi.org/10.1137/S0895480192243516

Motivation

David S. Johnson's 1974 paper Approximation Algorithms for Combinatorial Problems (J. Comput. System Sci. 9 (1974) 256–278) is one of the founding papers of the theory of approximation algorithms. It asks, for optimization problems whose decision versions Karp had just shown to be polynomial complete, how close a fast heuristic can be guaranteed to come to the optimum in the worst case, and it measures this with a worst-case performance ratio that is still the standard yardstick.

Its first example is SUBSET-SUM, the simplest form of the knapsack problem: pack items of given sizes into a knapsack of capacity $b$ so as to fill it as much as possible. For this problem the paper gives a family of algorithms $A_k$, one for each $k \ge 1$, whose guaranteed ratio $(k+1)/k$ tends to $1$. It is one of the first examples of what is now called a polynomial-time approximation scheme: for every $\epsilon > 0$ there is a polynomial-time algorithm within a factor $1 + \epsilon$ of optimal. Sahni (1975) extended the idea to the knapsack problem with utilities, and Ibarra and Kim (1975) later obtained fully polynomial schemes for knapsack and subset-sum.

This mission formalizes Theorem 1 of the paper, the performance guarantee of $A_k$ together with its tightness.

Setting

An input $\langle T, s, b\rangle$ of SUBSET-SUM is a finite set $T$, a positive rational size $s(x)$ for every $x \in T$, and a positive rational bound $b$. An approximate solution is a subset $T' \subseteq T$ with $m(T') \le b$, where the measure is $m(T') = \sum_{x \in T'} s(x)$. The problem is a maximization problem with optimal measure

$$\langle T, s, b\rangle^* = \max\{\, m(T') : T' \subseteq T,\ m(T') \le b \,\}.$$

Fix $k \ge 1$ and call $x$ big if $s(x) > b/(k+1)$ and small otherwise. Algorithm $A_k$ keeps a set $\mathrm{SUB}$, its measure $\mathrm{SUM}$, and the remaining elements $\mathrm{LEFT}$:

1. $\mathrm{SUB}$ is a subset of the big elements whose measure is as large as possible without exceeding $b$; $\mathrm{SUM} = m(\mathrm{SUB})$ and $\mathrm{LEFT} = T \setminus \mathrm{SUB}$.
2. If $s(x) + \mathrm{SUM} > b$ for every $x \in \mathrm{LEFT}$, return $\mathrm{SUB}$.
3. Otherwise pick $y \in \mathrm{LEFT}$ with $s(y) + \mathrm{SUM}$ as large as possible without exceeding $b$, move it from $\mathrm{LEFT}$ to $\mathrm{SUB}$, add $s(y)$ to $\mathrm{SUM}$, and return to step 2.

Steps 1 and 3 may have ties. Following the paper, a set $T_1$ is choosable by $A_k$ if some resolution of all ties produces it, and the performance $A_k(u)$ on input $u$ is the smallest measure of a choosable output. The ratio is $r(A_k, u) = u^*/A_k(u) \ge 1$, and $R[A_k](n)$ is its maximum over inputs of size at most $n$.

Formalization targets

Goal: Theorem 1 (p. 260)

For $k \ge 1$ and $n > 0$,

$$R[A_k](n) \le \frac{k+1}{k}, \qquad \lim_{n \to \infty} R[A_k](n) = \frac{k+1}{k}.$$

Formally, for every $k \ge 1$: every choosable output $T_1$ of every input satisfies $k\,\langle T,s,b\rangle^* \le (k+1)\,m(T_1)$; and for every $\delta > 0$ some input has a choosable output $T_1$ with $m(T_1) > 0$ and $\langle T,s,b\rangle^* > \big(\tfrac{k+1}{k} - \delta\big)\,m(T_1)$.

Milestones

1. For $T_1$ choosable and $T_0$ any approximate solution, $m(T_1^{\mathrm{BIG}}) \ge m(T_0^{\mathrm{BIG}})$ (p. 260).
2. If a small $x \in T$ is not in a choosable $T_1$, then $s(x) + m(T_1) > b$, hence $m(T_1) > kb/(k+1) \ge \tfrac{k}{k+1}\langle T,s,b\rangle^*$ (p. 261).
3. The stronger dichotomy: $m(T_1) = \langle T,s,b\rangle^*$ or $m(T_1) \ge \tfrac{k}{k+1}\,b$ (p. 260).
4. The lower-bound input $T = \{a_1,\dots,a_{k+2}\}$, $s(a_1) = 1+\varepsilon$, $s(a_i) = 1$ otherwise, $b = k+1$: its optimum is $k+1$, some output is choosable, and every choosable output has measure $k + \varepsilon$ (p. 261).

Significance

Theorem 1 shows that SUBSET-SUM admits polynomial-time algorithms with any worst-case ratio above $1$, in contrast with the other problems of the paper (set covering, graph colouring, maximum clique), whose best known ratios grow with the input. The algorithms $A_k$ are an early instance of the partial-enumeration schemes later used for knapsack-type problems. The tightness half shows that the analysis of $A_k$ itself cannot be sharpened.

The theorem has a short published proof, but no machine-checked version is known; there is no subset-sum or knapsack approximation result on the platform. The mission produces a reusable model of SUBSET-SUM, a model of nondeterministic algorithms through a run relation that captures every tie-break, and a checked proof that the worst case is exactly $(k+1)/k$. The same modelling pattern (choosable outputs, worst-case ratio taken over them) is used in the sibling missions of this series for MAX-SAT, set covering and exact covering.

Difficulty

The arithmetic of the upper bound is short; the difficulty is in reasoning about the algorithm as a nondeterministic process. The natural first attempt, implementing $A_k$ as a function with a fixed tie-breaking rule, proves a weaker statement: the guarantee must hold for every output the algorithm may return, including adversarial ties in step 1 (several maximum-measure sets of big elements) and step 3. Facts that are obvious for a single run, such as $\mathrm{SUM}$ always equalling $m(\mathrm{SUB})$ or which elements can enter $\mathrm{SUB}$ after step 1, have to be established for the run relation as a whole. The lower bound requires tracing the run on the explicit input for general $k$: exactly $k-1$ unit elements are added after $a_1$, and this must be shown for every choosable run, not only for one.

Formalization scope

• Numbers. Sizes and the bound are rationals (ℚ), as in the paper; sizes are required to be positive on $T$ and $b > 0$. The index $k$ is a natural number with $1 \le k$ as a hypothesis; $b/(k+1)$ is rational division, and "big" is the strict inequality $s(x) > b/(k+1)$.
• Optimum. opt u is Finset.sup' of the measure over the finite set of approximate solutions, which always contains $\emptyset$; it is $0$ when no element fits.
• Run relation. Choosable k u T₁ states that some admissible step 1 choice, followed by a finite chain of admissible iterations (Relation.ReflTransGen), reaches a halting state returning $T_1$. Every "closest to, without exceeding" is an existential choice among all maximizers.
• Size-free restatement. The paper's input size $|u|$ ("in some standard notation") is never fixed, so the goal quantifies over all inputs instead of over sizes. The upper bound for all choosable outputs is equivalent to $R[A_k](n) \le (k+1)/k$ for all $n$; since $R[A_k]$ is nondecreasing, the limit claim is equivalent to the supremum of the ratio over all inputs being $(k+1)/k$, which is the second part.
• Multiplicative ratios. No ratio is written as a division, so an output of measure $0$ cannot satisfy a bound vacuously; the lower-bound part requires $m(T_1) > 0$. The value $(k+1)/k$ is not claimed to be attained: the paper's family has ratio $(k+1)/(k+\varepsilon)$.
• Lower-bound input. A def on Fin (k + 2) exactly as on the page, with $0 < \varepsilon < 1$ (the page leaves the range implicit; $\varepsilon < 1$ keeps $a_1$ the only big element that fits when $k = 1$).
• Ruled out. A formalization with a deterministic tie-break, with a bound of the form $\mathrm{opt}/m \le c$ in a field where $x/0 = 0$, or with tightness for a single fixed $k$ would be trivial or weaker; none of these is the target.

Contributions welcome: proofs of the milestones and the goal, invariant lemmas for the run relation, and further sanity checks on small inputs. The running-time remark ($O(n^k)$ for step 1) and Sahni's knapsack extension are not part of the mission.

Selected references

• D. S. Johnson, Approximation algorithms for combinatorial problems, Journal of Computer and System Sciences 9 (1974) 256–278. https://doi.org/10.1016/S0022-0000(74)80044-9
• S. Sahni, Approximate algorithms for the 0/1 knapsack problem, Journal of the ACM 22 (1975) 115–124. https://doi.org/10.1145/321864.321873
• O. H. Ibarra, C. E. Kim, Fast approximation algorithms for the knapsack and sum of subset problems, Journal of the ACM 22 (1975) 463–468. https://doi.org/10.1145/321906.321909
• R. M. Karp, Reducibility among combinatorial problems, in Complexity of Computer Computations, Plenum (1972) 85–103. https://doi.org/10.1007/978-1-4684-2001-2_9

Motivation

A single machine must process a set of jobs, each with a processing time and a due-date, and a job that finishes after its due-date is late. Counting late jobs is the natural objective when a late order is simply lost, whatever its lateness. In the three-field notation of scheduling theory this is the problem $1\,\|\,\sum U_j$, and it is one of the few single-machine problems with a due-date objective that a simple greedy rule solves exactly.

J. Michael Moore gave that rule in 1968 (Management Science 15(1):102–109). The only exact method previously available was the Held–Karp dynamic program, which is exponential in the number of jobs. Moore's algorithm is two sorts plus at most $n(n+1)/2$ additions and comparisons. The rule, and the variant from the paper's Author's Supplement (credited to T. J. Hodgson and today called the Moore–Hodgson algorithm), is in every scheduling textbook, for example Brucker, Scheduling Algorithms, Ch. 4, and is the base case of later work on weighted and release-date variants.

Timeline:

• 1955: J. R. Jackson shows that a job set can be scheduled with no late job if and only if the earliest-due-date order has none (Management Science Research Project report 43, UCLA).
• 1968: Moore publishes the algorithm and its proof of optimality, with Hodgson's variant stated without proof.
• 1970s onward: the weighted version $1\,\|\,\sum w_jU_j$ is shown NP-hard (Karp 1972, via knapsack), and $1\,|\,r_j\,|\,\sum U_j$ likewise (Lenstra, Rinnooy Kan and Brucker 1977), so Moore's greedy rule does not extend to them.

Setting

A finite set $J$ of jobs is given. Job $j$ has a processing time $t_j \ge 0$ and a due-date $D_j$, and the paper assumes $t_j \le D_j$ for every job (a job that cannot finish on time even if started at time $0$ is removed beforehand). The machine starts at time $0$ and processes the jobs one after another, without idle time or preemption.

A schedule $S$ of $J$ is an ordering $(J_{i_1},\dots,J_{i_n})$ of all jobs of $J$. The job in position $k$ completes at $C_{i_k} = t_{i_1} + \dots + t_{i_k}$. The late set is $L = \{J_i : C_i > D_i\}$ and the early set is $E = \{J_i : C_i \le D_i\}$. A schedule is optimal if no schedule of $J$ has fewer late jobs. $A$ and $R$ denote the early and late jobs of $S$, each kept in their order in $S$.

Moore's algorithm works on a current sequence and a list of rejected jobs.

• Step 1: order the jobs by non-decreasing processing time (the shortest processing time rule).
• Step 2: find the first late job $J_{i_q}$ of the current sequence. If there is none, stop.
• Step 3: re-order $J_{i_1},\dots,J_{i_q}$ by non-decreasing due-date. If all of them are then early, keep the re-ordered sequence. Otherwise reject $J_{i_q}$ and remove it. Return to Step 2.

The output is the final current sequence sorted by due-dates, followed by the rejected jobs in any order.

In Lean, a schedule is IsSchedule J l, the late set is lateSet t D l, optimality is IsOptimal t D J l, $A$ and $R$ are earlyPart/latePart, and one pass of Steps 2–3 is the relation MooreStep t D, all in the namespace MooreLateJobs.NumLate.

Formalization targets

Goal: Moore's algorithm is optimal (The Algorithm, Step 2, p. 103)

Let $l_0$ be a shortest-processing-time schedule of $J$, and let a run of MooreStep from $(l_0,[\,])$ reach a state $(\mathrm{cur},\mathrm{rej})$ in which $\mathrm{cur}$ has no late job. Then for every due-date ordering $A_D$ of $\mathrm{cur}$ and every ordering $P$ of $\mathrm{rej}$,

$$(A_D,\,P)\ \text{is an optimal schedule for } J.$$

All tie-breaks in both sorts are covered.

Milestones

In attack order:

1. Lemma 1 (p. 105): every optimal schedule has the same number of late jobs as $(A,R)$ and as every $(A,P)$.
2. Jackson's lemma (p. 105).
3. Lemma 2 (p. 105): re-ordering $A$ by due-dates keeps an optimal $(A,R)$ schedule optimal.
4. Lemma 3 (p. 105): a job that is late in some optimal schedule can be removed and appended.
5. The repeated-elimination claim (p. 106): after removing jobs late in successive optimal schedules until the rest is feasible, $(A_D,P)$ is optimal.
6. Cases 2) and 3) of the Selection Algorithm (p. 107): in either case the job $J_q$ is late in some optimal schedule.
7. Progress and termination of the algorithm (p. 108).

A companion item states the p. 104 remark that the final current sequence need not be re-sorted: $(\mathrm{cur},P)$ is already optimal.

Significance

The theorem shows that the minimum number of late jobs on one machine can be found in $O(n\log n)$ time, by a rule that also produces an optimal schedule of a very particular shape: due-date ordered early jobs first, then the late jobs in any order. Lemma 3's decomposition, that jobs late in some optimal schedule may be discarded one at a time, is the template reused for many related greedy results in scheduling.

The result is classical and fully proved on paper. To our knowledge no machine-checked proof of Moore's algorithm, of the Moore–Hodgson variant, or of Jackson's rule exists in Mathlib. This mission produces a checked proof of the algorithm as stated in the paper, with every tie-break allowed, together with reusable single-machine objects (schedules as lists, completion times, late sets) and Jackson's earliest-due-date feasibility lemma.

Difficulty

Neither ordering rule works alone. Sorting by due-dates alone gives a schedule with no late job whenever one exists, but it can make many jobs late once any must be. Keeping the shortest jobs first does not respect the due-dates at all. The step that fails in a direct greedy argument is the claim that the specific job $J_{i_q}$, the one just found late, belongs to the late set of some optimal schedule. That job is not in general the longest job of the prefix, and the paper has to treat separately the two cases in which it is rejected. On top of this, the algorithm re-sorts prefixes on the fly, so the claim has to be tied to the invariants of the run: the prefix is early and due-date sorted, and the jobs after it are at least as long as $J_{i_q}$.

Formalization scope

• Jobs and times. Jobs form a type ι with decidable equality; $J$ is a Finset ι; $t, D : ι \to \mathbb{R}$.
• Standing hypotheses. Every statement that involves schedules assumes $t_j \ge 0$ and $t_j \le D_j$ on $J$. The first is added: processing times are durations, and Jackson's lemma fails for negative times. The second is the paper's assumption on p. 102.
• Schedules and completion times. A schedule is a duplicate-free list with exactly the jobs of $J$. Positions are 0-based, and the job in position $k$ completes at the sum of the first $k+1$ processing times. Lateness is strict ($C_j > D_j$).
• Optimality compares against every schedule of the same job set.
• Ties. Orderings "by due-dates" and "by processing times" are List.Pairwise with ≤. Ties are arbitrary, and every statement quantifies over all such orderings.
• The algorithm. Steps 2–3 are the relation MooreStep. The re-ordered prefix is any due-date sorted permutation of the first $q+1$ jobs, and case 2) rejects the first late job $J_{i_q}$ itself, not the longest job of the prefix (that is Hodgson's variant). A run is Relation.ReflTransGen.

The goal must concern runs of this step relation from a shortest-processing-time schedule of $J$. Replacing the run by an arbitrary set of rejected jobs satisfying invariants would state a different theorem. The goal is not vacuous: the progress and termination milestones show that a terminal state is always reached.

Contributions are welcome at every level. Useful ones include general lemmas on completion times under permutation and filtering of lists, a proof of Jackson's lemma, proofs of the Selection Algorithm cases, and a proof of Hodgson's variant.

Selected references

• J. M. Moore, An n Job, One Machine Sequencing Algorithm for Minimizing the Number of Late Jobs, Management Science 15(1):102–109, 1968. https://doi.org/10.1287/mnsc.15.1.102
• J. R. Jackson, Scheduling a Production Line to Minimize Maximum Tardiness, Research Report 43, Management Science Research Project, UCLA, 1955.
• M. Held and R. M. Karp, A Dynamic Programming Approach to Sequencing Problems, J. SIAM 10(1):196–210, 1962. https://doi.org/10.1137/0110015
• R. M. Karp, Reducibility among Combinatorial Problems, in Complexity of Computer Computations, 1972. https://doi.org/10.1007/978-1-4684-2001-2_9
• J. K. Lenstra, A. H. G. Rinnooy Kan and P. Brucker, Complexity of Machine Scheduling Problems, Annals of Discrete Mathematics 1:343–362, 1977. https://doi.org/10.1016/S0167-5060(08)70743-X
• P. Brucker, Scheduling Algorithms, 5th ed., Springer, 2007. https://doi.org/10.1007/978-3-540-69516-5

Motivation

A processor, a machine shop or a single server must work through a set of jobs one at a time, and each job is costly when it is late. Deciding the order is the single-machine sequencing problem, the simplest and most studied model of scheduling theory. Robert McNaughton's 1959 article Scheduling with Deadlines and Loss Functions (Management Science 6(1):1–12) treats it for a computer that must run several tasks, each with a deadline and a loss that grows linearly with the lateness. Its §2 gives the first sufficient condition under which a simple ratio rule is optimal in the presence of deadlines, and shows that interrupting and resuming tasks ("splitting", now called preemption) never helps on one processor.

Timeline.

• 1956: W. E. Smith, Various optimizers for single-stage production (Naval Research Logistics Quarterly 3), proves that sequencing jobs by non-increasing weight-to-processing-time ratio minimizes the total weighted completion time over non-preemptive sequences.
• 1959: McNaughton, §2 of the present paper, proves independently that the same ratio order is optimal against all schedules, split or not and with idle time (Theorem 2.3), and extends it to deadlines when no task finishes early in that order (Theorem 2.4). §3 of the same paper gives the "wrap-around" rule for preemptive makespan on identical processors, and §4 the non-preemptive optimality for weighted completion time on several processors.
• 1977: J. K. Lenstra, A. H. G. Rinnooy Kan and P. Brucker show that minimizing total weighted tardiness on one machine, the general problem of §2, is strongly NP-hard (Annals of Discrete Mathematics 1); this is why §2 gives a sufficient condition and not an algorithm.

Setting

There are $m$ tasks $(1),\dots,(m)$ for a single processor, and the present is time $0$. Task $(i)$ takes $a_i > 0$ units of processing time, has a deadline $d_i$ and a penalty rate $p_i \ge 0$. If $(i)$ is finished at time $C_i \le d_i$ there is no loss; otherwise the loss on $(i)$ is $p_i x$, where $x = C_i - d_i$ is the time from the deadline to the completion. Thus the loss on a task completed at time $t$ is

$$\ell_i(t) = p_i \max(0,\ t - d_i).$$

The ratio of task $(i)$ is $r_i = p_i / a_i$.

A task may be split: part of it may run between times 4 and 6 and the remainder between times 8 and 11, and similarly in any finite number of parts. A schedule $S$ is therefore a finite list of pieces, each a task together with a start and a stop time. It is feasible when every piece lies in $[0,\infty)$ with start $\le$ stop, no two pieces overlap in time, and the pieces of each task $(i)$ have total length exactly $a_i$. The completion time $C_i(S)$ is the latest stop time of a piece of $(i)$, and the total loss is

$$c(S) = \sum_{i=1}^{m} \ell_i\bigl(C_i(S)\bigr).$$

For an order $\sigma$ of the tasks ($\sigma(k)$ in position $k$), the sequenced schedule $S_\sigma$ runs the tasks without splits and without unused time: $\sigma(k)$ occupies $\bigl[\sum_{l<k} a_{\sigma(l)},\ \sum_{l\le k} a_{\sigma(l)}\bigr]$. The order is in decreasing $r_i$ when $k \le l$ implies $r_{\sigma(l)} \le r_{\sigma(k)}$. Finally $c^*(S)$ denotes the total loss of $S$ computed as if $d_1 = \dots = d_m = 0$.

Formalization targets

Goal: Theorem 2.4 (p. 5)

If $\sigma$ is in decreasing $r_i$ and no task finishes before its deadline in $S_\sigma$, i.e. $d_i \le C_i(S_\sigma)$ for every $i$, then $S_\sigma$ is feasible and

$$c(S_\sigma) \le c(S') \qquad \text{for every feasible schedule } S'.$$

The competitors $S'$ may split tasks and leave the processor idle. The condition is sufficient but not necessary.

Milestones, in attack order

1. Theorem 2.1 (p. 4): if both $(i)$ and $(j)$ run in the $a_i + a_j$ consecutive units of time after a time $t$ past both deadlines and $r_i > r_j$, their joint loss is strictly smaller when $(i)$ goes first:

$$\ell_i(t+a_i) + \ell_j(t+a_i+a_j) < \ell_j(t+a_j) + \ell_i(t+a_j+a_i).$$

2. The reduction in the proof of Theorem 2.2 (pp. 4–5): a feasible schedule with more than $m$ pieces can be replaced by a feasible one with fewer pieces and no greater loss.
3. Theorem 2.2 (p. 4): some optimal schedule, optimal among all feasible schedules, splits no task.
4. Theorem 2.3 (p. 5): if $d_1 = \dots = d_m = 0$, the sequenced schedule in decreasing $r_i$ minimizes the total loss over all feasible schedules.
5. The display of the proof of Theorem 2.4 (p. 6): if no task finishes early in $S = S_\sigma$, then for every feasible $S'$,

$$c(S') - c(S) \ge c^*(S') - c^*(S).$$

Significance

The result. Theorem 2.3 is the ratio rule for total weighted completion time, in its strongest single-machine form: it holds against preemptive schedules and schedules with idle time, not only against permutations. Theorem 2.4 carries the rule over to deadlines and linear tardiness penalties under a checkable condition on one schedule. Since weighted tardiness is strongly NP-hard in general, a condition of this kind is what one can hope for, and the paper's two-step heuristic for general deadlines (p. 6) is built on it. Theorem 2.2, as the paper remarks (p. 6), "does not depend on the linear loss function": it makes non-preemptive scheduling without loss of generality for single-machine objectives of this kind.

Formalizing it. All results of §2 are proved in the paper and are textbook material; none has a machine-checked proof on the platform. The platform's Scheduling Algorithms V mission formalizes the multi-processor results of §§3–4 (via Brucker's textbook), and nothing there states a single-processor ratio rule with deadlines. This mission supplies a single-processor schedule model with splitting, the interchange lemma, the non-preemption theorem and the ratio rule, each over all feasible schedules.

Difficulty

The interchange argument of Theorem 2.1 compares only two schedules that differ in the order of two adjacent tasks. Turning it into optimality against every feasible schedule requires two further steps, and each fails if done naively. First, a competitor may split tasks and leave gaps; the interchange argument does not apply to such schedules, so a separate argument must remove splits without raising any completion time. Second, with deadlines the loss $\max(0, t - d_i)$ is not linear in the completion time, so the ratio order is in general not optimal; the obvious attempt to repeat the interchange argument fails as soon as a task can finish before its deadline, since moving such a task later costs nothing. This is why Theorem 2.4 needs its hypothesis that no task finishes early, and why the paper leaves the general case to a heuristic.

Formalization scope

Tasks and positions are the zero-based indices of Fin m; times, lengths, deadlines and penalties are real numbers. A schedule is a List of pieces (task, start, stop), mirroring the public definition SchedulingAlgorithms_ParallelMachines with one processor. Feasibility requires $0 \le$ start $\le$ stop, pairwise disjoint pieces, and exact total length $a_i$ per task; zero-length pieces and unsorted lists are allowed. The completion time is the maximum stop time of the task's pieces ($0$ for a task with no pieces, which feasibility excludes). "No split" means exactly one piece per task, so two abutting pieces count as a split. "Decreasing $r_i$" is non-increasing, with ties in any order. "Minimal" and "optimal" are stated as $\le$ against every feasible schedule, never as an infimum.

Standing assumptions, stated in every item: $a_i > 0$ (tasks take time, and $r_i = p_i/a_i$ needs $a_i \ne 0$), and $p_i \ge 0$ for Theorems 2.2–2.4 and the proof steps (penalties are non-negative; with a negative penalty and idle time allowed the loss is unbounded below). Theorem 2.1 carries no sign condition. No condition is placed on the deadlines.

A formalization that restricts the competitors of Theorems 2.2–2.4 to unsplit schedules, or to sequenced schedules of other orders, states a weaker theorem and is ruled out: every statement quantifies over all feasible schedules.

A complete development needs: sums over sublists of pieces, rearrangements of pieces of a schedule and their effect on completion times, and optimality over permutations of a finite set of tasks. The schedule model and the non-preemption argument are reusable for any single-machine regular objective. Contributions of intermediate lemmas on these points are welcome.

Selected references

• R. McNaughton, Scheduling with Deadlines and Loss Functions, Management Science 6(1):1–12, 1959. https://doi.org/10.1287/mnsc.6.1.1
• W. E. Smith, Various optimizers for single-stage production, Naval Research Logistics Quarterly 3(1–2):59–66, 1956. https://doi.org/10.1002/nav.3800030106
• J. K. Lenstra, A. H. G. Rinnooy Kan, P. Brucker, Complexity of machine scheduling problems, Annals of Discrete Mathematics 1:343–362, 1977. https://doi.org/10.1016/S0167-5060(08)70743-X
• P. Brucker, Scheduling Algorithms, 5th ed., Springer, 2007. https://doi.org/10.1007/978-3-540-69516-5

Motivation

A semidefinite program (SDP) minimizes a linear objective $c^Tx$ subject to a linear matrix inequality $F(x) = F_0 + \sum_{i=1}^m x_i F_i \succeq 0$. In engineering applications the coefficient matrices are rarely known exactly: they come from measurements, from a model of a physical plant, or from a finite-precision implementation. El Ghaoui, Oustry and Lebret (SIAM J. Optim. 9(1), 1998) asked for solutions that remain feasible for every admissible value of the uncertain data, and showed how to compute such robust solutions by semidefinite programming. The paper appeared alongside Ben-Tal and Nemirovski's robust convex programming (Math. Oper. Res. 23(4), 1998) and is one of the two founding treatments of robust SDP.

When the uncertainty has structure (a block-diagonal perturbation, repeated scalar parameters, a symmetric matrix), the exact robust problem is NP-hard (El Ghaoui and Lebret, SIAM J. Matrix Anal. Appl. 18, 1997). This is the same obstacle that robust control meets in computing the structured singular value, and the remedy the paper uses, scaling matrices that commute with the perturbation structure, goes back to that literature (Doyle, IEE Proc. D 129, 1982; Fan, Tits and Doyle, IEEE Trans. Automat. Control 36, 1991). This mission formalizes the resulting tractable conservative approximation, Theorem 3.2 of the paper, together with the lemma it rests on and an application to integer feasibility problems.

Setting

Fix natural numbers $m, n, p, q$. The decision variable is $x \in \mathbb{R}^m$. The nominal data are affine maps

$$F(x) = F_0 + \sum_{i=1}^m x_i F_i \in \mathbb{R}^{n\times n}, \qquad R(x) = R_0 + \sum_{i=1}^m x_i R_i \in \mathbb{R}^{q\times n},$$

with every $F_i$ symmetric, and fixed matrices $L \in \mathbb{R}^{n\times p}$, $D \in \mathbb{R}^{q\times p}$. A perturbation is a matrix $\Delta \in \mathbb{R}^{p\times q}$, and the perturbed constraint matrix is the linear-fractional representation (LFR)

$$\mathbf{F}(x,\Delta) = F(x) + L\Delta(I - D\Delta)^{-1}R(x) + R(x)^T(I - \Delta^TD^T)^{-1}\Delta^TL^T,$$

which is defined when $\det(I - D\Delta) \neq 0$. The perturbation ranges over a linear subspace $\mathcal{D} \subseteq \mathbb{R}^{p\times q}$, which encodes the structure, and is bounded by a level $\rho > 0$ in the spectral norm $\|\Delta\|$ (the largest singular value). The robust feasible set is

$$\mathcal{X}_\rho = \{x : \text{for every } \Delta \in \mathcal{D} \text{ with } \|\Delta\| \le \rho,\ \det(I - D\Delta) \neq 0 \text{ and } \mathbf{F}(x,\Delta) \succeq 0\},$$

and the robust SDP (RSDP) is to minimize $c^Tx$ over $\mathcal{X}_\rho$.

The scaling set of $\mathcal{D}$ is the linear subspace

$$\mathcal{B} = \{(S,T,G) \in \mathbb{R}^{p\times p}\times\mathbb{R}^{q\times q}\times\mathbb{R}^{p\times q} : S\Delta = \Delta T,\ G\Delta^T = -\Delta G^T \text{ for every } \Delta \in \mathcal{D}\}.$$

Formalization targets

Goal: Theorem 3.2 (p. 37), as an inclusion of feasible sets

For every $x$: if some $(S,T,G) \in \mathcal{B}$ has $S \succ 0$, $T \succ 0$ and

$$\begin{bmatrix} F(x) - LSL^T & R(x)^T - LSD^T + LG \\ R(x) - DSL^T + G^TL^T & \rho^{-2}T - DSD^T + DG + G^TD^T\end{bmatrix} \succ 0,$$

then $x \in \mathcal{X}_\rho$, and in fact $\mathbf{F}(x,\Delta) \succ 0$ for every $\Delta \in \mathcal{D}$ with $\|\Delta\| \le \rho$. A companion item states the consequence for optimal values: the SDP value is an upper bound on the RSDP value, with both infima taken in the extended reals.

Milestones

1. Lemma 3.2 (p. 37): the same implication for constant $F$, $R$ and $\rho = 1$, with the matrix (13).
2. The full-perturbation case (p. 37): for $\mathcal{D} = \mathbb{R}^{p\times q}$ and $p, q \ge 1$, $\mathcal{B}$ consists exactly of the triples $(\tau I_p, \tau I_q, 0)$, with $\tau \ge 0$ when $S \succeq 0$.
3. Theorem 5.6 (p. 48): if $F_i = 2L_iR_i$ with $r_i = \operatorname{rank} F_i$, and $x_{\mathrm{feas}}$ satisfies, for some $\lambda \ge 0$ and block-diagonal $S = S^T$, $G = -G^T$,

$$\begin{bmatrix} F(x_{\mathrm{feas}}) - \lambda I - LSL^T & \tfrac12R^T + LG \\ \tfrac12R - GL^T & S\end{bmatrix} \succ 0,$$

then every integer vector closest to $x_{\mathrm{feas}}$ in the maximum norm satisfies $F(z) \succeq 0$.

Significance

The result. Theorem 3.2 replaces an NP-hard semi-infinite constraint, one matrix inequality for each admissible perturbation, by a single linear matrix inequality in the enlarged variable $(x, S, T, G)$. Every point it certifies is robustly feasible, so its optimal value is a certified upper bound on the robust optimum and its optimizer is a usable robust solution. In the full case the scalings collapse to one multiplier $\tau$ (milestone 2), which connects the bound to the exact reformulation of Section 3.1 of the paper. Theorem 5.6 shows the same machinery at work on a combinatorial problem: robustness against perturbations of size $1/2$ in each coordinate of $x$ turns an SDP-feasible point into an integer solution by rounding.

Formalizing it. The results are proved in the paper (Lemma 3.2 with the proof deferred to [16]); none of them has a machine-checked proof that this mission is aware of, and the platform has no linear-fractional or structured-perturbation results. The formalization also settles the exact form of the certificate: as printed, the matrix (13) and the LMI of Theorem 3.2 contain products that are dimensionally undefined, and this mission states the condition the proof actually yields (see the scope section).

Difficulty

The inequality to be proved is a statement about infinitely many perturbations, and $\mathbf{F}(x,\Delta)$ depends on $\Delta$ through a matrix inverse. The natural first step, eliminating $\Delta$ by an exact S-procedure as in the full case, is not available: with a structured $\mathcal{D}$ the set of pairs of vectors linked by some $\Delta \in \mathcal{D}$ is not described by one quadratic inequality, and losslessness fails. The scalings in $\mathcal{B}$ give several valid quadratic inequalities instead, and one must show that their combination controls every $\Delta$ in the norm ball, including the well-posedness claim $\det(I - D\Delta) \neq 0$, which is part of the conclusion rather than an assumption. The commutation condition $S\Delta = \Delta T$ must be turned into an inequality for $\|\Delta\| \le 1$, which requires more than the definition of the spectral norm. For Theorem 5.6 the block-diagonal perturbation family and the rescaling between $\rho = 1/2$ and the stated matrix must be matched to the general lemma.

Formalization scope

Matrices are Matrix (Fin a) (Fin b) ℝ; $\succ 0$ and $\succeq 0$ are Matrix.PosDef and Matrix.PosSemidef (both include symmetry); block matrices are Matrix.fromBlocks on Fin n ⊕ Fin q. The norm of a perturbation is the $\ell^2$ operator norm (open scoped Matrix.Norms.L2Operator), i.e. the largest singular value; the maximum norm in Theorem 5.6 is Mathlib's sup norm on Fin m → ℝ. $\mathcal{D}$ is a Submodule. Affine maps are given by coefficient families indexed by Fin (m+1). Mathlib's matrix inverse is $0$ at a singular matrix, so every statement pairs the LFR with $\det(I - D\Delta) \neq 0$. The standing assumption $\rho > 0$ of Section 3 is a hypothesis.

Readings and corrections of the printed statements:

• (13) as printed is dimensionally inconsistent; we state the condition the proof yields, which coincides with the printed one when $G$ is square and skew-symmetric and $\mathcal{D}$ consists of symmetric matrices. Concretely, (11) prints $G \in \mathbb{R}^{q\times p}$ with $G\Delta = -\Delta^TG^T$ and (13) prints the blocks $R - DSL - GL^T$ and $T - GD^T + DG - DSD^T$; the mission uses $G \in \mathbb{R}^{p\times q}$ with $G\Delta^T = -\Delta G^T$ and the blocks $R - DSL^T + G^TL^T$ and $T - DSD^T + DG + G^TD^T$. The same correction applies to the LMI of Theorem 3.2 (with $\rho^{-2}T$). Theorem 5.6 is stated as printed.
• "An upper bound on the RSDP (4) and a corresponding solution $x$ can be computed by solving the SDP" is read as the inclusion of the SDP's feasible projection in $\mathcal{X}_\rho$, for every $x$; the goal states it with the strict conclusion $\mathbf{F}(x,\Delta) \succ 0$ as well. The value form is a separate item.
• In the full-perturbation remark, "for some $\tau \ge 0$" is stated under $S \succeq 0$, and "We then recover the exact results of section 3.1" is not formalized.
• In Theorem 5.6, $\mathcal{S}$'s index range "$i = 1,\dots,n$" is read as $i = 1,\dots,m$; the hypothesis $r_i = \operatorname{rank}F_i$ is kept.

Trivializing formalizations are ruled out: $(0,0,0) \in \mathcal{B}$ always, so the hypotheses $S \succ 0$ and $T \succ 0$ are kept outside $\mathcal{B}$; $\mathcal{D}$ is a subspace, not an arbitrary set; and the norm is the spectral norm, not Mathlib's default entrywise norm.

A complete development needs the square root of a positive definite matrix and its commutation with $S$ and $T$, the spectral-norm characterization $\Delta\Delta^T \preceq \|\Delta\|^2 I$, Schur-complement and congruence facts for block matrices, and a linear-fractional identity relating $(I - D\Delta)^{-1}$ to an auxiliary vector. These are reusable well beyond this mission; contributions of any of them, and of the value and rounding corollaries, are welcome.

Selected references

• L. El Ghaoui, F. Oustry, H. Lebret, Robust Solutions to Uncertain Semidefinite Programs, SIAM J. Optim. 9(1):33–52, 1998. https://doi.org/10.1137/S1052623496305717
• L. El Ghaoui, H. Lebret, Robust solutions to least-squares problems with uncertain data, SIAM J. Matrix Anal. Appl. 18:1035–1064, 1997. https://doi.org/10.1137/S0895479896298130
• M. K. H. Fan, A. L. Tits, J. C. Doyle, Robustness in the presence of mixed parametric uncertainty and unmodeled dynamics, IEEE Trans. Automat. Control 36:25–38, 1991. https://doi.org/10.1109/9.62265
• J. C. Doyle, Analysis of feedback systems with structured uncertainties, IEE Proc. D 129(6):242–250, 1982. https://doi.org/10.1049/ip-d.1982.0053
• A. Ben-Tal, A. Nemirovski, Robust convex optimization, Math. Oper. Res. 23(4):769–805, 1998. https://doi.org/10.1287/moor.23.4.769
• S. Boyd, L. El Ghaoui, E. Feron, V. Balakrishnan, Linear Matrix Inequalities in System and Control Theory, SIAM, 1994. https://doi.org/10.1137/1.9781611970777

Motivation

Many firms learn about demand before it has to be served: customers place orders days or weeks ahead of the date they want delivery. Gallego and Özer (Management Science 47(10), 2001) model this advance demand information in a periodic-review inventory system and ask how the optimal replenishment policy should use it. Classical inventory theory (Arrow, Harris and Marschak 1951; Scarf 1959; Veinott 1965, 1966; Iglehart 1963) assumes that nothing about future demand is known when an order is placed. With advance orders the state of the system is no longer a single number, and it is not a priori clear whether the familiar policy structures survive.

This mission covers the paper's zero set-up cost case (Section 5). The companion mission Integrating Replenishment Decisions with Advance Demand Information I covers the positive set-up cost case and its $(s, S)$ policies.

Setting

Time is divided into periods $t = 1, \dots, T$. The supply lead time is an integer $L \ge 0$, and the information horizon is $N$. In period $t$ customers place orders $D_t = (D_{t,t}, \dots, D_{t,t+N})$, where $D_{t,s} \ge 0$ is demand placed in period $t$ for delivery in period $s$. Throughout, $N > L + 1$; write $M = N - L - 1 \ge 1$.

At the start of period $t$ the decision maker knows two things. The first is the modified inventory position $x_t$: on-hand stock plus outstanding orders minus backorders, net of the demand already observed for the protection period $t, \dots, t+L$. The second is the vector

$$o_t = (o_{t,t+L+1}, \dots, o_{t,t+N-1}) \in \mathbb{R}^M$$

of demands already observed for periods beyond the protection period. The decision maker raises the position to $y \ge x_t$ at zero fixed cost, the demand vector $D_t$ is realised, and the state moves to

$$x_{t+1} = y - \sum_{k=0}^{L+1} D_{t,t+k} - o_{t,t+L+1}, \qquad o_{t+1,s} = o_{t,s} + D_{t,s}\ (s = t+L+2, \dots, t+N),$$

with $o_{t,t+N} = 0$.

Costs enter through a single-period cost $G_t : \mathbb{R} \to \mathbb{R}$ (holding, backorder and linear ordering cost, charged against the demand over the protection period) and one-period discount factors $\alpha_{t+1} > 0$. The optimal cost-to-go $J_t$ and the cost $V_t$ of ordering up to $y$ satisfy

$$J_t(x, o) = \min_{y \ge x} V_t(y, o), \qquad V_t(y, o) = G_t(y) + \alpha_{t+1}\,\mathbb{E}\,J_{t+1}(x_{t+1}, o_{t+1}), \qquad J_{T+1} \equiv 0,$$

where the expectation is over $D_t$. The base-stock level in period $t$ is the smallest minimizer

$$y_t(o) = \min\{y : V_t(y, o) = \min_x V_t(x, o)\},$$

and the myopic level is the smallest minimizer of the single-period cost,

$$y^m_t = \min\{y : G_t(y) = \min_x G_t(x)\}.$$

A function $f(x, \theta)$ has decreasing differences if $f(x_1, \theta) - f(x_2, \theta) \le f(x_1, \theta') - f(x_2, \theta')$ whenever $x_1 \ge x_2$ and $\theta \ge \theta'$ componentwise.

Formalization targets

Goal: Theorem 5 (p. 1352)

If $t \mapsto y^m_t$ is nondecreasing on $\{1, \dots, T\}$, then for every period $t$ and every observed-demand vector $o \ge 0$,

$$y_t(o) = y^m_t .$$

The optimal order-up-to level then ignores all advance information beyond the protection period. A second item states the paper's stationary special case: if $G_t = G$ for all $t$, the smallest minimizer $y^m$ of $G$ is the optimal base-stock level in every period.

Milestones: Theorem 4 (p. 1351)

For every period $t$ and every fixed $o_t$:

1. $V_t(\cdot, o_t)$ is convex and $V_t(x, o_t) \to \infty$ as $|x| \to \infty$;
2. $y_t(o_t)$ exists and $J_t(x, o_t) = V_t(\max(y_t(o_t), x), o_t)$: a state-dependent base-stock policy is optimal;
3. $J_t(\cdot, o_t)$ is nondecreasing and convex;
4. $V_t(x, o)$ has decreasing differences in $(x, o)$;
5. $J_t(x, o)$ has decreasing differences in $(x, o)$;
6. $y_t(o)$ is nondecreasing in $o$.

Parts 1–3 are what the goal's proof uses. Parts 4–6 are the paper's second zero set-up result, monotonicity of the base-stock level in observed demand.

Significance

The theorem identifies when advance demand information beyond the protection period can be ignored. When the myopic levels do not decrease over time, which includes stationary costs and ramping-up demand, the $(1 + M)$-dimensional dynamic program collapses to a sequence of one-dimensional newsvendor-type problems. That is both a computational simplification and a managerial statement: information about demand after the protection period does not change the order. Theorem 4, Part 5 gives the complementary monotone comparative statics. When the myopic condition fails, more observed demand never lowers the order-up-to level.

The results are proved in the paper, with the proofs in Appendix B. No machine-checked version exists. This mission produces a formal account of the finite-horizon recursion with a multi-dimensional information state, and checks the base-stock and myopic-optimality arguments against it.

Difficulty

The obvious induction carries convexity of $J_{t+1}$ backward, but here the future cost is evaluated at a random next state $(x_{t+1}, o_{t+1})$ whose first coordinate depends on the current observed demand $o_{t,t+L+1}$. Showing that the base-stock level does not depend on $o_t$ therefore needs more than convexity. It needs to know where $J_{t+1}(\cdot, o_{t+1})$ is flat, uniformly in the random $o_{t+1}$, and that the next position cannot exceed the current order-up-to level. The latter holds only on the reachable states, where observed demands are nonnegative. For a sufficiently negative $o_{t,t+L+1}$ the next period starts above its myopic level whatever is ordered now, and the conclusion fails. On the analytic side, every infimum and expectation in the recursion must be shown to be finite and attained before the order-theoretic argument can start.

Formalization scope

The model is parametrised by $L$ and $M \ge 1$, with $N = L + M + 1$. The demand vector is a function on $\{0, \dots, N\}$ and $o$ a function on $\{0, \dots, M-1\}$, ordered componentwise. $J_t$ is defined by backward recursion with $J_t \equiv 0$ for $t > T$. The minimum over $y \ge x$ is a real infimum and the expectation a Bochner integral against the law $\mu_t$ of $D_t$. Attainment and finiteness are consequences proved in the theorems, not assumptions. Base-stock and myopic levels are characterised as smallest minimizers (IsLeast), never through sInf.

The single-period cost $G_t$ is a primitive rather than being assembled from $c_t$, $g_t$ and the lead-time demand; the paper's $G_t$ has the assumed properties, so the theorems cover the paper's model. Hypotheses the paper uses without stating, all placed on primitives and labelled in the statements:

• nonnegative demands, $D_{t,s} \ge 0$ almost surely;
• coercivity of $G_t$ (the paper states it for $\tilde G_t$ only);
• $\alpha_{t+1} > 0$;
• finiteness of the expectation in (9), guaranteed by linear growth of $G_t$ and finite first moments of $D_t$. This covers piecewise-linear holding and backorder costs with any finite-mean demand (including the paper's Poisson example), but excludes superlinear costs;
• in the goal, $o \ge 0$, the set of reachable states.

The goal cannot be trivialised: the hypotheses are satisfied by concrete instances (for example $G_t(y) = |y|$ with any finite-mean nonnegative demand), and the conclusion identifies the base-stock level exactly rather than asserting that some minimizer exists.

Out of scope: the reduction of the control problem to the functional equation (Appendix A, Özer 2000), the infinite-horizon Theorem 6, and Lemma 5, whose proof argues on the integers and whose real-valued form with a unit forward difference is unverified. Contributions of general lemmas are welcome: convexity and attainment for $\inf_{y \ge x}$ of a convex coercive function, and preservation of convexity and decreasing differences under expectation. All of them are reusable in other inventory models.

Selected references

• G. Gallego, Ö. Özer, Integrating Replenishment Decisions with Advance Demand Information, Management Science 47(10):1344–1360, 2001. https://doi.org/10.1287/mnsc.47.10.1344.10261
• A. F. Veinott, Optimal Policy for a Multi-Product, Dynamic, Nonstationary Inventory Problem, Management Science 12(3):206–222, 1965. https://doi.org/10.1287/mnsc.12.3.206
• D. L. Iglehart, Optimality of (s, S) Policies in the Infinite Horizon Dynamic Inventory Problem, Management Science 9(2):259–267, 1963. https://doi.org/10.1287/mnsc.9.2.259
• D. M. Topkis, Supermodularity and Complementarity, Princeton University Press, 1998. https://doi.org/10.1515/9781400822539